System and method for generating and protecting cryptographic keys

ABSTRACT

In the present disclosure, implementations of Diffie-Hellman key agreement are provided that, when embodied in software, resist extraction of cryptographically sensitive parameters during software execution by white-box attackers. Four embodiments are taught that make extraction of sensitive parameters difficult during the generation of the public key and the computation of the shared secret. The embodiments utilize transformed random numbers in the derivation of the public key and shared secret. The traditional attack model for Diffie-Hellman implementations considers only black-box attacks, where attackers analyze only the inputs and outputs of the implementation. In contrast, white-box attacks describe a much more powerful type of attacker who has total visibility into the software implementation as it is being executed.

RELATED APPLICATION DATA

This application claims priority to International Patent Application No. PCT/CA2012/000114, filed Feb. 9, 2012, the disclosure of which is hereby incorporated by reference in its entirety

FIELD OF THE INVENTION

The present invention relates to methods of generating and protecting cryptographic keys, and is particularly concerned with doing so in a white-box attack environment.

BACKGROUND OF THE INVENTION

Traditionally, cryptographic implementations were designed to resist only black-box attacks. These are attacks where an adversary attempts to deduce the cryptographic key inside the implementation by analyzing only inputs and outputs. In a black-box attack, it is assumed that the adversary knows what family of cryptographic algorithm they are targeting (e.g., AES, DES, RSA, etc.), but all other details (e.g. execution time, power consumption, memory accesses) are unavailable to them. In essence, a black-box attacker treats the implementation as a mathematical function. However, a white-box attacker is a much more powerful type of adversary and can analyze all parts of the implementation. Rather than just study inputs and outputs, a white-box attacker can see everything that goes on inside the implementation. For example, if the attacker is targeting cryptographic software running on, say, a PC or mobile phone, then they can execute that software inside a debugger and examine memory and register values during its execution. Conventional software implementations of cryptographic algorithms are completely insecure against white-box attackers since the cryptographic key can be observed in memory. New implementations of cryptographic algorithms that resist white-box attacks have recently been proposed (e.g., AES [1], DES [2], ECDSA [3], SHA [4]); however, there have not yet been any proposals for protecting key agreement algorithms, such as Diffie-Hellman [5].

The systems and methods disclosed herein provide techniques for implementing a key agreement scheme that resists white-box attacks.

SUMMARY OF THE INVENTION

An object of the present invention is to provide methods for protecting cryptographic keys and other sensitive parameters from a white-box attacker in implementations of key agreement schemes such as Diffie-Hellman.

Elliptic curve Diffie-Hellman (ECDH) is a variant of traditional Diffie-Hellman (DH) where computations are done in an algebraic group defined by points on an elliptic curve. ECDH is defined in the document “SEC 1: Elliptic Curve Cryptography, Version 2.0” [8] and in “Suite B Implementer's Guide to NIST SP 800-56A” [7]. Introductory information on elliptic curve cryptography and ECDH can be found in the text “Guide to Elliptic Curve Cryptography” by Hankerson, Menezes and Vanstone [9]. At high level, ECDH works as follows: two parties exchange public keys and combine them with their respective private keys to derive a shared secret. Even though their respective private keys are different, the mathematics underlying ECDH ensures that each party will derive the same shared secret. Private keys are typically just random numbers, and a corresponding public key is computed from a private key using a mathematical operation that is computationally easy to evaluate but very computationally expensive to invert.

ECDH is utilized in a number of cryptographic protocols including Transport Layer Security (TLS), which is used to secure Internet web browsing, and the Full Authentication protocol, which is part of the Digital Transmission Content Protection (DTCP) system [10]. The shared key established using ECDH can be used to build a secure channel between end-parties (i.e. using symmetric key cryptography for message confidentiality and authentication). Any attacker sitting between the end-parties that records and analyzes messages exchanged during the ECDH scheme should not be able to recover the shared key unless they carry out some computation which is widely believed to be infeasible. However, in some applications, the attacker does not sit between the end-parties; instead, the attacker is able to observe the ECDH computations taking place at one of the end-parties. This is a potential concern in DTCP.

DTCP is a digital rights management (DRM) scheme and is commonly used to enforce permissions on video content in home networks. For example, a user might wish to stream video from their set-top box to a personal device like a laptop or mobile phone. If the rights associated with the video content permit this activity, the set-top box will initiate ECDH with the personal device to derive a shared key; the shared key is used to encrypt video content as it is transmitted to the personal device where it will be played. If an attacker was able to determine the shared key, then it would be able to record the video stream and decrypt it outside the DRM system, thereby freeing it from any restrictions. In this scenario, a legitimate user may act as an attacker if they are motivated to free video content from the DRM system. Users are able to observe the computations that take place on their personal device using a debugger and other tools. For this reason, malicious users can be modeled as white-box attackers. This leads to the following question: can an ECDH shared secret be computed in a white-box environment without exposing the shared secret to an attacker?

In the present disclosure, four embodiments of ECDH that are resistant to white-box attacks are provided; these embodiments can be used to build software implementations that are able to derive cryptographic keys and establish trust in hostile execution environments. Although our description focuses on the elliptic curve variant of Diffie-Hellman, one skilled in the art will appreciate that these embodiments can easily be adapted to any Diffie-Hellman scheme; this includes, in particular, Diffie-Hellman schemes where the underlying algebraic group is a subgroup of a finite field.

The four embodiments generate and manipulate transformed random numbers. A transformed random number is an obfuscated or masked form of an original random number. In each embodiment, a transformed random number, denoted by {circumflex over (r)} or {circumflex over ({circumflex over (r)}, is used to generate an ECDH public key. The same transformed random number is then used to generate the shared secret. The resultant public key, generated using the transformed random number, will be the same as what would be generated using the original random number. This property also holds for generation of the shared secret. To accomplished this, some auxillary elliptic curve points must be pre-computed and utilized in the computations.

In the first embodiment, we generate a transformed random number {circumflex over (r)} using a linear transformation defined by two instance-specific, compile time, random numbers k₁ and k₂ chosen from the interval 1 to n−1 where n denotes the size of the underlying elliptic curve group. In the second embodiment, we generate a transformed random number {circumflex over (r)} using a linear transformation defined by one instance-specific, compile time, random number k₁, chosen from the interval 1 to n−1. This approach is similar to the first embodiment except that only one compile time random number is used to generate the transformed random number {circumflex over (r)}. In the third embodiment, a transformed random number {circumflex over ({circumflex over (r)} is generated by applying a pseudo-Hadamard transformation [11] to the transformed random number {circumflex over (r)} that is generated in the first embodiment. In the fourth embodiment, a transformed random number {circumflex over ({circumflex over (r)} is generated applying a pseudo-Hadamard transformation to the transformed-random number {circumflex over (r)} that is generated in the second embodiment.

BRIEF DESCRIPTION OF THE DRAWINGS

The present invention will be further understood from the following detailed description with reference to the drawings in which:

FIG. 1 illustrates public key generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with a first embodiment of the present disclosure;

FIG. 2 illustrates shared secret generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with the first embodiment of the present disclosure;

FIG. 3 illustrates public key generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with a second embodiment of the present disclosure;

FIG. 4 illustrates shared secret generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with the second embodiment of the present disclosure;

FIG. 5 illustrates public key generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with a third embodiment of the present disclosure;

FIG. 6 illustrates shared secret generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with the third embodiment of the present disclosure;

FIG. 7 illustrates public key generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with a fourth embodiment of the present disclosure; and

FIG. 8 illustrates shared secret generation for White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) in accordance with the fourth embodiment of the present disclosure.

SYMBOLS USED TO DESCRIBE THE PREFERRED EMBODIMENT

The following symbols are used within this disclosure to describe embodiments of the present disclosure.

Symbols Description q An ECC domain parameter; the field size. FR Field Representation indicator. An indication of the basis used for representing field elements. For the Suite B curves, FR is NULL. a, b ECC domain parameters; two field elements that define the equation of an elliptic curve. SEED An ECC domain parameter; an initialization value that is used during domain parameter generation that can also be used to provide assurance at a later time that the resulting domain parameters were generated arbitrarily. G An ECC domain parameter, which is a distinguished point on an elliptic curve that generates the subgroup of order n. h An ECC domain parameter; the cofactor, which is the order of the elliptic curve divided by the order of the point G. For the Suite B curves, h = 1. n An ECC domain parameter; the order of the point G. D The set of ECC domain parameters, (q, FR, a, b{, SEED}, G, n, h). d_(e, A), d_(e, B) Party A's and Party B's ephemeral private keys. These are integers in the range [1, n − 1]. d_(s, A), d_(s, B) Party A's and Party B's static private keys. These are integers in the range [1, n − 1]. O The point at infinity; a special point in an elliptic curve group that serves as the (additive) identity. Q_(e, A), Q_(e, B) Party A's and Party B's ephemeral public keys. These are points on the elliptic curve defined by the domain parameters. Q_(s, A), Q_(s, B) Party A's and Party B's static public keys. These are points on the elliptic curve defined by the domain parameters. x_(p), y_(p) Elements of the finite field of size q, representing, respectively, the x and y coordinates of a point P. For Suite B curves, these are integers in the interval [0, q − 1]. Z A shared secret that is used to derive secret keying material using a key derivation function.

Elliptic Curve Diffie-Hellman Schemes

ECDH is recommended by the NSA in their Suite B cryptography guidelines [6]. More precisely, Suite B recommends the use of Ephemeral Unified ECDH (EU-ECDH) and One-Pass ECDH (OP-ECDH). These two ECDH schemes are described in the Suite B implementer's guide [7]. In the Ephemeral Unified scheme, each party generates an ephemeral key pair each time they want to compute a new shared secret. In the One-Pass scheme, an ephemeral key pair is generated by only one party; the other party uses a static key pair, which may be reused in the computation of many shared secrets. The implementer's guide states that EU-ECDH is preferred over OP-ECDH; however, there are some scenarios where EU-ECDH cannot be used. This is the case, for example, in store-and-forward communication systems where one party is not available to contribute an ephemeral public key. In scenarios like this, OP-ECDH can be used.

Exposure of Sensitive Parameters

In both the EU-ECDH and OP-ECDH schemes, there are two main phases of computation: key pair generation and shared secret computation. The steps required in each phase are described in detail in the implementer's guide [7]; however, the computations in those steps leak sensitive parameters to white-box attackers.

The procedure for key pair generation in [7] specifies that the key pair (private key d_(A) and public key Q_(A)) be generated based on a random number r as follows: d_(A)=r+1, and Q_(A)=d_(A)G. Note that “d_(A)G” denotes the elliptic curve point formed by adding the generator G to itself d_(A) times. The procedure for shared secret computation specifies that the shared secret, Z, be computed as the x-coordinate of an elliptic curve point P defined as follows: P=d_(A)Q_(B)=(r+1)Q_(B), where d_(A) is party A's private key and Q_(B) is party B's public key.

Straightforward implementations of key pair generation and shared secret computation on party A's device will expose the secret key d_(A) in that device's memory. This is apparent for the computations that produce d_(A)=r+1, Q_(A)=d_(A)G, and P=d_(A)Q_(B). Thus, a white-box attacker that examines memory-dumps of party A's device will be able to deduce the private key, d_(A), as well as the shared secret, Z. To maintain security against white-box attackers, the computations executed for key pair generation and shared secret computation must be protected.

DETAILED DESCRIPTION OF THE PREFERRED EMBODIMENT

FIG. 1 illustrates the process of key pair generation for our proposed White-Box Elliptic Curve Diffie-Hellman (WB-ECDH) implementation, in accordance with a first embodiment of the present disclosure. FIG. 2 illustrates the process of shared secret computation.

Key Pair Generation.

Each static and ephemeral private key d and public key Q shall be generated using the appropriate domain parameters, D, which must be selected from one of the two sets of domain parameters specified in the list of Suite B elliptic curves and domain parameters (see Appendix A of [7]). Henceforth, we assume that domain parameters have been appropriately selected.

Target:

Compute Q=dG without disclosing d inside the memory of one party's machine, where Q=dG=(r+1)G, r is the random number. Q is the public key, and may represent any one of Q_(e,A), Q_(e,B) or Q_(s,B) depending on the particular ECDH scheme (EU-ECDH or OP-ECDH). d is the private key, and, similarly, it may represent any one of d_(e,A), d_(e,B) or d_(s,B).

Transformed Random Number:

In this embodiment, the transformation t of a random number r is denoted as {circumflex over (r)} and defined as {circumflex over (r)}=t(r)=k₁r+k₂ mod n, where k₁ and k₂ are two instance-specific, compile time, random numbers and k₁, k₂εZ_(n)* (note that Z_(n)* denotes the set of integers in the range 1 to n−1 that are relatively prime to the domain parameter n). Thus, {circumflex over (r)}=k₁r+k₂ mod n and r=k₁ ⁻¹{circumflex over (r)}−k₂k₁ ⁻¹ mod n. It is important to note that the output of the random bit generator (RBG) used in key pair generation is interpreted as producing transformed random numbers. This means that the transformation t(r) does not have to be explicitly applied by the implementation—the transformation is implicit in the way in which the output of the RBG is interpreted.

Precomputation:

Once the instance-specific, compile time, random numbers k₁ and k₂ are selected, two auxiliary elliptic curve points are pre-computed: Ĝ=k₁ ⁻¹G and U₂=(k₁−k₂)Ĝ. These two points are computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time).

The steps of key pair generation are as follows:

Input:

None.

Output:

-   -   1. Status: The status returned from the key pair generation         procedure. The status will indicate SUCCESS or an ERROR.     -   2. {circumflex over (r)} A transformed random number.     -   3. Q: The generated public key.

Process:

-   -   1. Set N=len(n) (i.e. set N equal to the bit length of the         domain parameter n). Check that

N is valid, that is, N=256 or N=384 (the only valid lengths for Suite B).

-   -   2. If N is invalid, then return an ERROR indication.     -   3. Set requested_security_strength=the security strength         associated with N (either 128 when using P-256 or 192 when using         P-384).     -   4. Obtain a string of N bits (b′₁, b′₂, . . . , b′_(N)) from a         Random Bit Generator (RBG) with a security strength of         requested_security_strength or more, which generates the         transformed random number {circumflex over (r)}=k₁r+k₂ mod n. If         an ERROR indication is returned, then return the ERROR         indication.     -   5. Compute U₁={circumflex over (r)}Ĝ (recall that Ĝ=k₁ ⁻¹G was         precomputed).

6. Compute Q=U₁+U₂ (recall that U₂=(k₁−k₂)Ĝ was precomputed).

-   -   7. Return SUCCESS, {circumflex over (r)} and Q.

Verification:

It can be shown that the previous procedure produces a public key that is compatible with the original procedure given in [7]. The original procedure produces a public key Q=(r+1)G; for the new procedure, we have

$\begin{matrix} {Q = {U_{1} + U_{2}}} \\ {= {{\hat{r}\hat{G}} + {\left( {k_{1} - k_{2}} \right)\hat{G}}}} \\ {= {{\hat{r}k^{- 1}G} + {\left( {k_{1} - k_{2}} \right)k_{1}^{- 1}G}}} \\ {= {\left( {{k_{1}^{- 1}\hat{r}} - {k_{2}k_{1}^{- 1}} + {k_{1}k_{1}^{- 1}}} \right)G}} \\ {= {\left( {{k_{1}^{- 1}\hat{r}} - {k_{2}k_{1}^{- 1}} + 1} \right)G}} \\ {= {\left( {r + 1} \right)G}} \end{matrix}$

Cost:

In our proposed key pair generation procedure, the cost of elliptic curve computations is as follows. Three elliptic curve point multiplications are required, two of which are offline and one is online. One elliptic curve point addition is required.

Shared Secret Computation.

Party A computes the shared secret Z, which is the x-coordinate of an elliptic curve point, using the appropriate domain parameters D, party B's public key (Q_(B)), and their transformed random number {circumflex over (r)}_(A). Party B computes the shared secret in an analogously, using party A's public key and their own transformed random number. We continue

Target:

Compute P=(r_(A)+1)Q_(B) without disclosing r_(A) (party A's random number) inside the memory of party A's machine.

Precomputation:

Let k_(1(A)) and k_(2(A)) denote party A's instance-specific, compile time, random numbers. Party A should pre-compute the integer values s=k_(1(A)) ⁻¹ mod n and t=k_(1(A))−k_(2(A)) mod n. These two values are computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time). If party B's public key is known in advance, as is the case in OP-ECDH, then further pre-computation can be done. When doing OP-ECDH, the elliptic curve points {circumflex over (Q)}_(B) ₁ =sQ_(B)=k_(1(A)) ⁻¹Q_(B) and {circumflex over (Q)}_(B) ₃ =t{circumflex over (Q)}_(B) ₁ =(k_(1(A))−k_(2(A))){circumflex over (Q)}_(B) ₁ should be computed off-line.

The steps of shared secret computation are as follows:

Input:

-   -   1. {circumflex over (r)}_(A): Party A's transformed random         number (recall that r_(A)=k_(1(A)) ⁻¹{circumflex over         (r)}_(A)−k_(2(A))k_(1(A)) ⁻¹ mod n).     -   2. Q_(B): The other party's (party B) public key.

Output:

The shared secret Z or an error indicator.

Process:

-   -   1. Compute points {circumflex over (Q)}_(B) ₁ =sQ_(B)=k_(1(A))         ⁻¹Q_(B); {circumflex over (Q)}_(B) ₂ ={circumflex over         (r)}_(A){circumflex over (Q)}_(B) ₁ ; and {circumflex over         (Q)}_(B) ₃ =t{circumflex over (Q)}_(B) ₁         =(k_(1(A))−k_(2(A))){circumflex over (Q)}_(B) ₁     -   2. Compute P={circumflex over (Q)}_(B) ₂ +{circumflex over         (Q)}_(B) ₃     -   3. If P=O, the point at infinity, output an error indicator.     -   4. Z=x_(P) where x_(P) is the x-coordinate of P.

Verification:

It can be shown that the previous procedure, which is summarized in FIG. 2, produces a shared secret that is compatible with the original procedure given in [7]. The original procedure produces a shared secret equal to (r_(A)+1)Q_(B); for the new procedure, we have

$\begin{matrix} {P = {{\hat{Q}}_{B_{2}} + {\hat{Q}}_{B_{3}}}} \\ {= {{{\hat{r}}_{A}{\hat{Q}}_{B_{1}}} + {\left( {k_{1{(A)}} - k_{2{(A)}}} \right){\hat{Q}}_{B_{1}}}}} \\ {= {{{\hat{r}}_{A}k_{1{(A)}}^{- 1}Q_{B}} - {k_{2{(A)}}k_{1{(A)}}^{- 1}Q_{B}} + {k_{1{(A)}}k_{1{(A)}}^{- 1}Q_{B}}}} \\ {= {\left( {{k_{1{(A)}}^{- 1}{\hat{r}}_{A}} - {k_{2{(A)}}k_{1{(A)}}^{- 1}} + {k_{1{(A)}}k_{1{(A)}}^{- 1}}} \right)Q_{B}}} \\ {= {\left( {r_{A} + 1} \right)Q_{B}}} \end{matrix}$

Cost:

In our proposed shared secret computation procedure, the cost of elliptic curve computations is as follows. Three elliptic curve point multiplications required. For OP-ECDH, two of the point multiplications are off-line and one is online; for EU-ECDH, all three point multiplications are online. One elliptic curve point addition is required.

The second embodiment of the disclosure is a reduced version of the first embodiment. The second embodiment is described with reference to FIGS. 3 and 4. The main concern in this design is computational efficiency.

Key Pair Generation.

We now describe key pair generation.

Target:

Compute Q=dG without disclosing d inside the memory of one party's machine.

Transformed Random Number:

In this embodiment, the transformation of the random number r is defined as {circumflex over (r)}=t(r)=k₁r mod n, where k₁ is an instance specific, compile time random number selected from the set Z_(n)*. Thus, {circumflex over (r)}=k₁r mode and r=k₁ ⁻¹{circumflex over (r)} mod n.

Precomputation:

Once the instance-specific, compile time, random number k₁ is selected, an auxiliary elliptic curve point is pre-computed: Ĝ=k₁ ⁻¹G. As in the first embodiment, this point should be computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time).

The steps of key pair generation are as follows:

Input and Output are the same as that of the first embodiment.

Process:

-   -   1. Steps 1, 2 and 3 are the same as of the first embodiment.     -   4. Obtain a string of N bits (b′₁, b′₂, . . . , b′_(N)) from a         Random Bit Generator (RBG) with a security strength of         requested_security_strength or more, which generates the         transformed random number {circumflex over (r)}=k₁r mod n. If an         ERROR indication is returned, then return the ERROR indication.     -   5. Compute U₁={circumflex over (r)}Ĝ (recall that Ĝ=k₁ ⁻¹G was         pre-computed).     -   6. Compute Q=U₁+G     -   7. Return SUCCESS, {circumflex over (r)} and Q.

Verification:

The following sequence of identities show that the previous procedure produces a public key that is compatible with the original procedure given in [7]:

$\begin{matrix} {Q = {U_{1} + G}} \\ {= {{\hat{r}*\hat{G}} + G}} \\ {= {{\hat{r}*k_{1}^{- 1}*G} + G}} \\ {= {\left( {{k_{1}^{- 1}*\hat{r}} + 1} \right)*G}} \\ {= {\left( {r + 1} \right)G}} \end{matrix}$

Cost:

The costs of elliptic curve computations for key pair generation are now as follows. Two elliptic curve point multiplications are required (one online and one offline). One elliptic curve point addition is required.

Shared Secret Computation:

We explain shared secret computation from the perspective of Party A. Party A will use the appropriate domain parameters D, the other party's public key (Q_(B)), and their transformed random number {circumflex over (r)}_(A) to compute Z, which is the x-coordinate of an elliptic curve point P.

Target:

Compute P=(r_(A)+1)Q_(B) without disclosing r_(A) inside the memory of party A's machine.

Precomputation:

Let k_(1(A)) denote party A's instance-specific, compile time, random number. Party A should pre-compute the integer values s=k_(1(A)) ⁻¹ mod n. This value should be computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time). If party B's public key is known in advance, as is the case in OP-ECDH, then further pre-computation can be done. When doing OP-ECDH, the elliptic curve point {circumflex over (Q)}_(B)=sQ_(B)=k_(1(A)) ⁻¹Q_(B) should be computed off-line.

The steps are as follows:

Input:

-   -   1. {circumflex over (r)}_(A): Party A's transformed random         number (recall that r_(A)=k_(1(A)) ⁻¹{circumflex over (r)}_(A)         mod n).     -   2. Q_(B): The other party's (party B) public key.

Output:

The shared secret Z or an error indicator.

Process:

-   -   1. Compute points {circumflex over (Q)}_(B) ₁ =sQ_(B)=k_(1(A))         ⁻¹Q_(B); and {circumflex over (Q)}_(B) ₂ ={circumflex over         (r)}_(A){circumflex over (Q)}_(B) ₁     -   2. P={circumflex over (Q)}_(B) ₂ +Q_(B)     -   3. If P=O, the point at infinity, output an error indicator.     -   4. Z=x_(P) where x_(P) is the x-coordinate of P.

Verification:

It can be shown that the previous procedure, which is summarized in FIG. 4, produces a shared secret that is compatible with the original procedure given in [7]. Consider the following sequence of identities:

$\begin{matrix} {P = {{\hat{Q}}_{B_{2}} - Q_{B}}} \\ {= {{{\hat{r}}_{A}*{\hat{Q}}_{B_{1}}} + Q_{B}}} \\ {= {{{\hat{r}}_{A}*k_{1{(A)}}^{- 1}*Q_{B}} + Q_{B}}} \\ {= {\left( {{k_{1{(A)}}^{- 1}*{\hat{r}}_{A}} + 1} \right)*Q_{B}}} \\ {= {\left( {r_{A} + 1} \right)*Q_{B}}} \end{matrix}$

Cost:

The costs of elliptic curve computations for shared secret computation are now as follows. Two elliptic curve point multiplications are required. For OP-ECDH, one point multiplication is off-line and one is online. For EU-ECDH, both point multiplications are online. One elliptic curve point addition is required (for both OP-ECDH and EU-ECDH).

The second embodiment is a reduced version of the first embodiment. One instance specific random number is removed from the first embodiment, which reduces the number of elliptic curve computations, but possibly sacrifices security strength. This methodology is suitable for those applications where computational efficiency is a prime concern.

The third embodiment is described with reference to FIGS. 5 and 6.

Key Pair Generation:

The third embodiment is an enhanced version of the first embodiment where security strength is a prime concern. The computational cost of this embodiment is higher.

Target:

Compute Q=dG without disclosing d inside the memory of one party's machine.

Transformed Random Numbers:

In this methodology, the Random Bit Generator generates two transformed random numbers {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ that, in combination, encode two random numbers r_(i) and r₂. The resultant public key, Q, may be generated using either r₁ or r₂, and this choice may be made at runtime within the key pair generation procedure. Thus, our proposed procedure can compute Q=dG=(r₁+1)G or Q=dG=(r₂+1)G. The values {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ output by the RBG are interpreted as being produced by two layers of transformations. One layer corresponds to the same transformation technique of the first embodiment; namely, a linear transform using instance specific, compile time random numbers drawn from the set Z_(n)*. The second layer is a pseudo-Hadamard transformation [11], which ensures that r₁ and r₂ can only be recovered from the combination of both {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂.

More precisely, {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ are interpreted as follows:

{circumflex over ({circumflex over (r)}₁₁ ={circumflex over (r)} ₁ +{circumflex over (r)} ₂ mod 2^(N)

{circumflex over ({circumflex over (r)}₂₂ ={circumflex over (r)} ₁+2{circumflex over (r)} ₂ mod 2^(N)

We interpret {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ as being the result of a pseudo-Hadamard transform applied to the values {circumflex over (r)}₁ and {circumflex over (r)}₂. Note that N is equal to the bit length of the domain parameter n, and {circumflex over (r)}₁ and {circumflex over (r)}₂ are transformed values that encode r₁ and r₂ respectively. To recover {circumflex over (r)}₂ from {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂, we compute {circumflex over ({circumflex over (r)}₂₂−{circumflex over ({circumflex over (r)}₁₁ mod 2^(N); and to recover {circumflex over (r)}₁ we compute 2{circumflex over ({circumflex over (r)}₁₁−{circumflex over ({circumflex over (r)}₂₂ mod 2^(N). The definition of {circumflex over (r)}₁ and {circumflex over (r)}₂ is

{circumflex over (r)} ₁ =k ₁₁ r ₁ +k ₁₂ mod n

{circumflex over (r)} ₂ =k ₂₁ r ₂ +k ₂₂ mod n

Here, k₁₁, k₁₂, k₂₁, k₂₂ are instance specific, compile time random numbers drawn from the set Z_(n)*. As in the first embodiment, we have r₁=k₁₁ ⁻¹{circumflex over (r)}₁−k₁₂k₁₁ ⁻¹ mod n and r₂=k₂₁ ⁻¹{circumflex over (r)}₂−k₂₂k₂₁ ⁻¹ mod n.

Precomputation:

Once the instance-specific, compile time, random numbers k₁₁, k₁₂, k₂₁, k₂₂ are selected, two sets of auxiliary elliptic curve points are pre-computed: Ĝ₁₁=k₁₁ ⁻¹G, Ĝ₁₂=k₁₁ ⁻¹Ĝ₁₁, U₁₂=(k₁₁−k₁₂)Ĝ₁₁ and Ĝ₂₁=k₂₁ ⁻¹G, Ĝ₂₂=k₂₁ ⁻¹Ĝ₂₁, U₂₂=(k₂₁−k₂₂)Ĝ₂₁. As in the first embodiment, these points should be computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time).

The steps of key pair generation are as follows:

Input and Output are the same as that of the first embodiment.

Process:

-   -   1. Steps 1, 2 and 3 are the same as that of the first         embodiment.     -   4. Obtain strings of N bits (b′₁₁, b′₁₂, . . . , b′_(IN)) and         (b′₂₁, b′₂₂, . . . , b′_(2N)) from a Random Bit Generator (RBG)         with a security strength of requested_security_strength or more,         which generates the transformed random numbers {circumflex over         ({circumflex over (r)}₁₁=({circumflex over (r)}₁+{circumflex         over (r)}₂)mod 2^(N) and {circumflex over ({circumflex over         (r)}₂₂=({circumflex over (r)}₁+2{circumflex over (r)}₂)mod         2^(N). If an ERROR indication is returned, then return the ERROR         indication.     -   5. If (“r₁” is chosen), then         -   a. Compute u₁₁=(2{circumflex over ({circumflex over             (r)}₁₁−{circumflex over ({circumflex over (r)}₂₂ mod             2^(N))k₁₁ mod n.         -   b. Compute Q=u₁₁Ĝ₁₂+U₁₂ (recall that Ĝ₁₂ and U₁₂ were             pre-computed)     -   6. If (“r₂” is chosen), then         -   c. Compute u₂₁=({circumflex over ({circumflex over             (r)}₂₂−{circumflex over ({circumflex over (r)}₁₁             mod^(2N))k₂₁.         -   d. Q=u₂₁*Ĝ₂₂+U₂₂ (recall that Ĝ₂₂ and U₂₂ were pre-computed)     -   7. Return SUCCESS, {circumflex over ({circumflex over (r)}₁₁,         {circumflex over ({circumflex over (r)}₂₂ and Q.

Verification:

The procedure is illustrated in FIG. 5 and it can be verified that the resultant Q is equal to either (r₁+1)G or (r₂+1)G as follows:

$\begin{matrix} {Q = {{u_{11}{\hat{G}}_{12}} + U_{12}}} \\ {= {{\left( {{2\; {\hat{\hat{r}}}_{11}} - {{\hat{\hat{r}}}_{22}{mod}\; 2^{N}}} \right)k_{11}k_{11}^{- 1}*{\hat{G}}_{11}} + U_{12}}} \\ {= {{{\hat{r}}_{1}k_{11}^{- 1}G} + {\left( {k_{11} - k_{12}} \right){\hat{G}}_{11}}}} \\ {= {\left( {{k_{11}^{- 1}{\hat{r}}_{1}} - {k_{12}k_{11}^{- 1}} + {k_{11}k_{11}^{- 1}}} \right)G}} \\ {= {\left( {r_{1} + 1} \right)G}} \end{matrix}$ $\begin{matrix} {Q = {{u_{21}*{\hat{G}}_{22}} + U_{22}}} \\ {= {{\left( {{\hat{\hat{r}}}_{22} - {{\hat{\hat{r}}}_{11}{mod}\; 2^{N}}} \right)*k_{21}*k_{21}^{- 1}*{\hat{G}}_{21}} + U_{22}}} \\ {= {{{\hat{r}}_{2}*k_{21}^{- 1}*G} + {\left( {k_{21} - k_{22}} \right)*{\hat{G}}_{21}}}} \\ {= {\left( {{k_{21}^{- 1}*{\hat{r}}_{2}} - {k_{22}*k_{21}^{- 1}} + {k_{21}*k_{21}^{- 1}}} \right)*G}} \\ {= {\left( {r_{2} + 1} \right)G}} \end{matrix}$

Shared Secret Computation:

We explain shared secret computation from the perspective of Party A. Party A will use the appropriate domain parameters D, the other party's public key (Q_(B)), and their own transformed random numbers {circumflex over ({circumflex over (r)}_(11(A)),{circumflex over ({circumflex over (r)}_(22(A)) to compute Z.

Target:

Without disclosing r_(1(A)) and r_(2(A)) inside the memory of the party A's machine, compute P=(r_(1(A))+1)Q_(B) or P=(r_(2(A))+1)Q_(B) depending on whether r_(1(A))+1 or r_(2(A))+1 was selected as party A's private key.

Precomputation:

Let k_(11(A)), k_(12(A)), k_(21(A)), k_(22(A)) denote party A's instance-specific, compile time, random numbers. Party A should pre-compute the integer values s₁=k_(11(A)) ⁻¹ mod n and s₂=k_(21(A)) ⁻¹ mod n. These values should be computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time). If party B's public key is known in advance, as is the case in OP-ECDH, then further pre-computation can be done. When doing OP-ECDH, the elliptic curve points {circumflex over (Q)}_(B) ₁₁ =s₁Q_(B)=k_(11(A)) ⁻¹Q_(B), {circumflex over (Q)}_(B) ₁₂ =(k_(11(A))−k_(12(A))){circumflex over (Q)}_(B) ₁₁ m {circumflex over (Q)}_(B) ₂₁ =s₂Q_(B)=k_(21(A)) ⁻¹Q_(B), and {circumflex over (Q)}_(B) ₂₂ =(k_(21(A))−k_(22(A))){circumflex over (Q)}_(B) ₂₁ should be computed off-line.

The steps are as follows:

Input:

-   -   1. {circumflex over ({circumflex over (r)}_(11(A)),{circumflex         over ({circumflex over (r)}_(22(A)): Party A's transformed         random numbers.     -   2. Q_(B): The other party's (party B) public key.

Output:

The shared secret Z or an error indicator.

Process:

-   -   1. If (“r₁” was chosen), then         -   A. Compute {circumflex over (Q)}_(B) ₁₁ =s₁Q_(B),             {circumflex over (Q)}_(B) ₁₂             =(k_(11(A))−k_(12(A))){circumflex over (Q)}_(B) ₁₁ and             {circumflex over (Q)}_(B) ₁₃ =(2{circumflex over             ({circumflex over (r)}_(11(A))−{circumflex over ({circumflex             over (r)}_(22(A))mod 2^(N)){circumflex over (Q)}_(B) ₁₁             (recall that s₁ was pre-computed).         -   B. P={circumflex over (Q)}_(B) ₁₂ +{circumflex over (Q)}_(B)             ₁₃         -   C. If P=O, the point at infinity, output an error indicator.         -   D. Z=x_(P) where x_(P) is the x-coordinate of P.     -   2. If (“r₂” was chosen), then         -   A. Compute {circumflex over (Q)}_(B) ₂₁ =s₂Q_(B),             {circumflex over (Q)}_(B) ₂₂             =(k_(21(A))−k_(22(A))){circumflex over (Q)}_(B) ₂₁ and             {circumflex over (Q)}_(B) ₂₃ =({circumflex over ({circumflex             over (r)}_(22(A))−{circumflex over ({circumflex over             (r)}_(11(A))mod 2^(N)){circumflex over (Q)}_(B) ₂₁ (recall             that s₂ was pre-computed).         -   B. P={circumflex over (Q)}_(B) ₂₂ +{circumflex over (Q)}_(B)             ₂₃         -   C. If P=O, the point at infinity, output an error indicator.         -   D. Z=x_(P) where x_(P) is the x-coordinate of P.

Verification:

The procedure is illustrated in FIG. 6. The correctness of the procedure can be verified as follows:

$\begin{matrix} {P = {{\hat{Q}}_{B_{12}} + {\hat{Q}}_{B_{13}}}} \\ {= {{\hat{Q}}_{B_{12}} + {\left( {{2*{\hat{\hat{r}}}_{11{(A)}}} - {{\hat{\hat{r}}}_{22{(A)}}{mod}\mspace{14mu} 2^{N}}} \right)*{\hat{Q}}_{B_{11}}}}} \\ {= {\left( {{k_{11{(A)}}^{- 1}*{\hat{r}}_{1{(A)}}} - {k_{12{(A)}}*k_{11{(A)}}^{- 1}} + {k_{11}*k_{11}^{- 1}}} \right)*Q_{B}}} \\ {= {\left( {r_{1{(A)}} + 1} \right)Q_{B}}} \end{matrix}$ $\begin{matrix} {P = {{\hat{Q}}_{B_{22}} + {\hat{Q}}_{B_{23}}}} \\ {= {{\hat{Q}}_{B_{22}} + {\left( {{\hat{\hat{r}}}_{22{(A)}} - {{\hat{\hat{r}}}_{11{(A)}}{mod}\mspace{14mu} 2^{N}}} \right)*{\hat{Q}}_{B_{21}}}}} \\ {= {\left( {{k_{21{(A)}}^{- 1}*{\hat{r}}_{2{(A)}}} - {k_{22{(A)}}*k_{21{(A)}}^{- 1}} + {k_{21}*k_{21}^{- 1}}} \right)*Q_{B}}} \\ {= {\left( {r_{2{(A)}} + 1} \right)Q_{B}}} \end{matrix}$

The third embodiment may be the most resistant to white-box attacker. In this embodiment, two random numbers use to make it more difficult for the attacker to follow the execution path of the program. Furthermore, two transformations are used to mask the random numbers. As a result, computational efficiency is degraded. The third embodiment is suitable for those applications where security strength is the main concern.

The fourth embodiment is described with reference to FIGS. 7 and 8.

Key Pair Generation:

The fourth embodiment is an enhanced version of the second embodiment.

Target:

Compute Q=dG without disclosing d inside the memory of one party's machine.

Transformed Random Numbers:

As was done the in third embodiment, the fourth embodiment uses the random bit generator to generate two transformed random numbers {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ that, in combination, encode two random numbers r₁ and r₂. The resultant public key, Q, may be generated using either r₁ or r₂, and this choice may be made at runtime within the key pair generation procedure. The values {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ output by the RBG are interpreted as being produced by two layers of transformations. One layer corresponds to the transformation technique of the second embodiment; namely, a linear transform using instance specific, compile time random numbers drawn from the set Z_(n)*. The second layer is a pseudo-Hadamard transformation, as in the third embodiment.

More precisely, {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ are interpreted as follows:

{circumflex over ({circumflex over (r)}₁₁ ={circumflex over (r)} ₁ +{circumflex over (r)} ₂ mod 2^(N)

{circumflex over ({circumflex over (r)}₂₂ ={circumflex over (r)} ₁+2{circumflex over (r)} ₂ mod 2^(N)

We interpret {circumflex over ({circumflex over (r)}₁₁ and {circumflex over ({circumflex over (r)}₂₂ as being the result of a pseudo-Hadamard transform applied to the values {circumflex over (r)}₁ and {circumflex over (r)}₂. Note that N is equal to the bit length of the domain parameter n, and {circumflex over (r)}₁ and {circumflex over (r)}₂ are transformed values that encode r₁ and r₂ respectively. To recover {circumflex over (r)}₂ from {circumflex over (r)}₁₁ and {circumflex over (r)}₂₂, we compute {circumflex over ({circumflex over (r)}₂₂−{circumflex over ({circumflex over (r)}₁₁ mod 2^(N); and to recover {circumflex over (r)}₁ we compute 2{circumflex over ({circumflex over (r)}₁₁−{circumflex over ({circumflex over (r)}₂₂ mod 2^(N). The definition of {circumflex over (r)}₁ and {circumflex over (r)}₂ is

{circumflex over (r)} ₁ =k ₁₁ r ₁ mod n

{circumflex over (r)} ₂ =k ₂₁ r ₂ mod n

Here, k₁₁, k₂₁ are instance specific, compile time random numbers drawn from the set Z_(n)*. As in the second embodiment, we have r₁=k₁₁ ⁻¹{circumflex over (r)}₁ mod n and r₂=k₂₁ ⁻¹{circumflex over (r)}₂ mod n.

Precomputation:

Once the instance-specific, compile time, random numbers k₁₁,k₂₁ are selected, two sets of auxiliary elliptic curve points are pre-computed: Ĝ₁₁=k₁₁ ⁻¹G, Ĝ₁₂=k₁₁ ⁻¹Ĝ₁₁ and Ĝ₂₁=k₂₁ ⁻¹G, Ĝ₂₂=k₂₁ ⁻¹Ĝ₂₁. As in the first embodiment, these points should be computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time).

The steps of key pair generation are as follows:

Input and Output are the same as that of the third embodiment.

Process:

-   -   1. Steps 1, 2 and 3 are the same as that of the first         embodiment.     -   4. Obtain strings of N bits (b′₁₁, b′₁₂, . . . , b′_(1N)) and         (b′₂₁, b′₂₂, . . . , b′_(2N)) from a Random Bit Generator (RBG)         with a security strength of requested_security_strength or more,         which generates the transformed random numbers {circumflex over         ({circumflex over (r)}₁₁=({circumflex over (r)}₁+{circumflex         over (r)}₂)mod 2^(N) and {circumflex over ({circumflex over         (r)}₂₂=({circumflex over (r)}₁+2{circumflex over (r)}₂)mod         2^(N). If an ERROR indication is returned, then return the ERROR         indication, and Invalid_Q.     -   5. If (“r₁” is chosen), then         -   a. Compute u₁₁=(2{circumflex over ({circumflex over             (r)}₁₁−{circumflex over ({circumflex over (r)}₂₂ mod             2^(N))k₁₁         -   b. Compute Q=u₁₁Ĝ₁₂+G (recall that Ĝ₁₂ was pre-computed).     -   6. If (“r₂” is chosen), then         -   a. Compute u_(2i)=({circumflex over ({circumflex over             (r)}₂₂−{circumflex over ({circumflex over (r)}₁₁ mod             2^(N))k₂₁         -   b. Q=u₂₁*Ĝ₂₂+G (recall that Ĝ₂₂ was pre-computed).     -   7. Return SUCCESS, {circumflex over ({circumflex over (r)}₁₁,         {circumflex over ({circumflex over (r)}₂₂ and Q.

Verification:

The computational process is illustrated in FIG. 7. Its correctness can be verified as follows:

$\begin{matrix} {Q = {{u_{11}*{\hat{G}}_{12}} + G}} \\ {= {{\left( {{2*\; {\hat{\hat{r}}}_{11}} - {{\hat{\hat{r}}}_{22}{mod}\; 2^{N}}} \right)*k_{11}*k_{11}^{- 1}*{\hat{G}}_{11}} + G}} \\ {= {{{\hat{r}}_{1}*k_{11}^{- 1}*G} + G}} \\ {= {\left( {{k_{11}^{- 1}*{\hat{r}}_{1}} + 1} \right)*G}} \\ {= {\left( {r_{1} + 1} \right)G}} \end{matrix}$ $\begin{matrix} {Q = {{u_{21}*{\hat{G}}_{22}} + G}} \\ {= {{\left( {{\hat{\hat{r}}}_{22} - {{\hat{\hat{r}}}_{11}{mod}\; 2^{N}}} \right)*k_{21}*k_{21}^{- 1}*{\hat{G}}_{21}} + G}} \\ {= {{{\hat{r}}_{2}*k_{21}^{- 1}*G} + G}} \\ {= {\left( {{k_{21}^{- 1}*{\hat{r}}_{2}} + 1} \right)*G}} \\ {= {\left( {r_{2} + 1} \right)G}} \end{matrix}$

Shared Secret Computation.

As before, we explain shared secret computation from the perspective of Party A. Party A will use the appropriate domain parameters D, the other party's public key (Q_(B)), and their own transformed random numbers {circumflex over ({circumflex over (r)}_(11(A)),{circumflex over ({circumflex over (r)}_(22(A)) to compute Z.

Target:

Without disclosing r_(1(A)) and r_(2(A)) inside the memory of the party A's machine, compute P=(r_(1(A))+1)Q_(B) or P=(r_(2(A))+1)Q_(B) depending on whether r_(1(A))+1 or r_(2(A))+1 was selected as party A's private key.

Precomputation:

Let k_(11(A)), k_(21(A)) denote party A's instance-specific, compile time, random numbers. Party A should pre-compute the integer values s₁=k_(11(A)) ⁻¹ mod n and s₂=k_(21(A)) ⁻¹ mod n. These values should be computed during an off-line set-up phase where the implementation is not subject to white-box attacks (e.g. at compilation time). If party B's public key is known in advance, as is the case in OP-ECDH, then further pre-computation can be done. When doing OP-ECDH, the elliptic curve points {circumflex over (Q)}_(B) ₁₁ =s₁Q_(B)=k_(11(A)) ⁻¹Q_(B) and {circumflex over (Q)}_(B) ₂₁ =s₂Q_(B)=k_(21(A)) ⁻¹Q_(B) should be computed off-line.

The steps of shared secret computation are as follows:

Input:

-   -   1. {circumflex over ({circumflex over (r)}_(11(A)),{circumflex         over ({circumflex over (r)}_(22(A)): Party A's transformed         random numbers.     -   2. Q_(B): The other party's (party B) public key.

Output:

The shared secret Z or an error indicator.

Process:

-   -   1. If (“r₁” was chosen), then         -   A. Compute {circumflex over (Q)}_(B) ₁₁ =s₁Q_(B) and             {circumflex over (Q)}_(B) ₁₂ =(2{circumflex over             ({circumflex over (r)}_(11(A))−{circumflex over ({circumflex             over (r)}_(22(A))mod 2^(N)){circumflex over (Q)}_(B) ₁₁         -   B. Compute P={circumflex over (Q)}_(B) ₁₂ +Q_(B)         -   C. If P=O, the point at infinity, output an error indicator.         -   D. Z=x_(P) where x_(P) is the x-coordinate of P.     -   2. If (“r₂” was chosen), then         -   A. Compute {circumflex over (Q)}_(B) ₂₁ =s₂Q_(B); and             {circumflex over (Q)}_(B) ₂₂ =({circumflex over ({circumflex             over (r)}_(22(A))−{circumflex over ({circumflex over             (r)}_(11(A)) mod 2^(N)){circumflex over (Q)}_(B) ₂₁         -   B. Compute P={circumflex over (Q)}_(B) ₂ +Q_(B)         -   C. If P=O, the point at infinity, output an error indicator.         -   D. Z=x_(P) where x_(P) is the x-coordinate of P.

Verification:

The computational process is illustrated in FIG. 8. Its correctness can be verified as follows:

$\begin{matrix} {P = {{\hat{Q}}_{B_{12}} + Q_{B}}} \\ {= {{\left( {{2*{\hat{\hat{r}}}_{11{(A)}}} - {{\hat{\hat{r}}}_{22{(A)}}{mod}\mspace{14mu} 2^{N}}} \right)*{\hat{Q}}_{B_{11}}} + Q_{B}}} \\ {= {\left( {{k_{11{(A)}}^{- 1}*{\hat{r}}_{1{(A)}}} + 1} \right)*Q_{B}}} \\ {= {\left( {r_{1{(A)}} - 1} \right)Q_{B}}} \end{matrix}$ $\begin{matrix} {P = {{\hat{Q}}_{B_{22}} + Q_{B}}} \\ {= {{\left( {{\hat{\hat{r}}}_{22{(A)}} - {{\hat{\hat{r}}}_{11{(A)}}{mod}\mspace{14mu} 2^{N}}} \right)*{\hat{Q}}_{B_{21}}} + Q_{B}}} \\ {= {\left( {{k_{21{(A)}}^{- 1}*{\hat{r}}_{2{(A)}}} + 1} \right)*Q_{B}}} \\ {= {\left( {r_{2{(A)}} + 1} \right)Q_{B}}} \end{matrix}$

The fourth embodiment attempts to reach a compromise between security and computational efficiency.

Table A shows the computational comparison among the proposed White-Box ECDH design embodiments. Table A allows the four embodiments to be ranked in terms of their computational cost. The embodiments can also be approximately ranked in descending order of security strength as embodiment 3, embodiment 1, embodiment 4, embodiment 2.

Numerous modifications, variations and adaptations may be made to the particular embodiments described above without departing from the scope patent disclosure, which is defined in the claims.

TABLE A Computational comparison among the proposed White-Box ECDH design methodologies Multiplication Addition Subtraction (Elliptic Curve Point) (Elliptic Curve Point) (Scalar) KEY PAIR EMBODIMENT 1 1 (online) 1 (online) 1 GENERATION 2 (offline) EMBODIMENT 2 1 (online) 1 (online) 0 1 (offline) EMBODIMENT 3 1 (online) 1 (online) 2 3 (offline) Scalar: 2 (for r₁) and 1 (for r₂) EMBODIMENT 4 1 (online) 1 (online) 1 2 (offline) Scalar: 2 (for r₁) and 1 (for r₂) STANDARD BLACK-BOX 1 0 0 SHARED SECRET EMBODIMENT 1 OP- ECDH 1 (online) 1 (online) 1 COMPUTATION 2 (offline) EU- ECDH 3 (online) 1 (online) 1 EMBODIMENT 2 OP- ECDH 2 (offline) 1 (online) 0 EU- ECDH 2 (online) 1 (online) 0 EMBODIMENT 3 OP- ECDH 1 (online) 1 (online) 2 2 (offline) Scalar: 1 (for r₁) and 0 (for r₂) EU- ECDH 3 (online) 1 (online) 2 Scalar: 1 (for r₁) and 0 (for r₂) EMBODIMENT 4 OP- ECDH 1 (online) 1 (online) 1 1 (offline) Scalar: 1 (for r₁) and 0 (for r₂) EU- ECDH 2 (online) 1 (online) 1 Scalar: 1 (for r₁) and 0 (for r₂) STANDARD BLACK- OP- ECDH 1 (offline) 0 0 BOX EU- ECDH 1 (online) 0 0

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1. A method of generating public key for an elliptic curve Diffie-Hellman (ECDH) key exchange protocol without disclosing the private key during computation of the public key, the method comprising the steps of: generating a random number; interpreting that random number as a linearly transformed random number; and generating a public key using the transformed number.
 2. The method of claim 1, wherein the linear transformation is {circumflex over (r)}=k₁r+k₂ mod n where k₁ and k₂ are two instance specific, compile time random values and are elements of Z*_(n).
 3. The method of claim 2 wherein the step of generating a public key computes Ĝ=k₁ ⁻¹G, U₁={circumflex over (r)}Ĝ, and U₂=(k₁−k₂)Ĝ and then Q=U₁+U₂.
 4. The method of claim 2 wherein k₂=0.
 5. The method of claim 4 wherein the step of generating a public key computes Ĝ=k₁ ⁻¹G, U₁={circumflex over (r)}Ĝ, and then Q=U₁+G.
 6. The method of claim 1 wherein the step of generating a random number generates two random numbers r₁ and r₂.
 7. The method of claim 6 wherein the step of linearly transforming the random number includes the steps of first applying a linear transform to random numbers r₁ and r₂, then applying a second transformation thereto.
 8. The method of claim 7 wherein the linear transformation is, r ₁=k₁₁r₁+k₁₂ mod n and {circumflex over (r)}₂=k₂₁r₂+k₂₂ mod n, where k₁₁, k₁₂, k₂₁, k₂₂ are instance specific, compile time randomly picked numbers and k₁₁, k₁₂, k₂₁, k₂₂εZ_(n)*.
 9. The method of claim 8 wherein the second transformation is t₂ on the transformed-random numbers {circumflex over (r)}₁ and {circumflex over (r)}₂ and is defined as {circumflex over ({circumflex over (r)}₁₁=({circumflex over (r)}₁+{circumflex over (r)}₂)mod 2^(N), and {circumflex over ({circumflex over (r)}₂₂=({circumflex over (r)}₁+2{circumflex over (r)}₂)mod 2^(N).
 10. The method of claim 9 wherein is r₁ chosen for computing the public key.
 11. The method of claim 10 where the step of generating a public key computes Ĝ₁₁=k₁₁ ⁻¹Ĝ; Ĝ₁₂=k₁₁ ⁻¹Ĝ₁₁; u₁₁=(2{circumflex over ({circumflex over (r)}₁₁−{circumflex over ({circumflex over (r)}₂₂ mod 2^(N))k₁₁ mod n and U₁₂=(k₁₁−k₁₂)Ĝ₁₁ and Q=u₁₁Ĝ₁₂+U₁₂.
 12. The method of claim 9 wherein is r₂ chosen for computing the public key.
 13. The method of claim 12 where the step of generating a public key computes Ĝ₂₁=k₂₁ ⁻¹G, Ĝ₂₂=k₂₁ ⁻¹Ĝ₂₁, u₂₁=({circumflex over ({circumflex over (r)}₂₂−{circumflex over ({circumflex over (r)}₁₁ mod 2^(N))k₂₁ mod n and U₂₂=(k₂₁−k₂₂)Ĝ₂₁ and Q=u₂₁*Ĝ₂₂+U₂₂.
 14. The method of claim 7 wherein the linear transformation is, {circumflex over (r)}₁=k₁₁r₁ and {circumflex over (r)}=k₂₁r₂, where k₁₁, and k₂₁ are instance specific, compile time randomly picked numbers, and k₁₁,k₂₁εZ_(n)*. Thus, r₁=k₁₁ ⁻¹{circumflex over (r)}₁, and r₂=k₂₁ ⁻¹{circumflex over (r)}₂.
 15. The method of claim 14 wherein the second transformation on the transformed-random numbers {circumflex over (r)}₁ and {circumflex over (r)}₂ is defined as {circumflex over ({circumflex over (r)}₁₁=({circumflex over (r)}₁+{circumflex over (r)}₂)mod 2^(N), and {circumflex over ({circumflex over (r)}=({circumflex over (r)}₁+2{circumflex over (r)}₂)mod 2^(N).
 16. The method of claim 15 wherein is r₁ chosen for computing the public key.
 17. The method of claim 16 wherein the step of generating a public key computes Ĝ₁₁=k₁₁ ⁻¹G; Ĝ₁₂=k₁₁ ⁻¹Ĝ₁₁; u₁₁=(2{circumflex over ({circumflex over (r)}₁₁−{circumflex over ({circumflex over (r)}₂₂ mod 2^(N))k₁₁ mod n; and Q=u₁₁Ĝ₁₂+G.
 18. The method of claim 15 wherein is r₂ chosen for computing the public key.
 19. The method of claim 18 wherein the step of generating a public key computes Ĝ₂₁=k₂₁ ⁻¹G; Ĝ₂₂=k₂₁ ⁻¹Ĝ₂₁; u₂₁=({circumflex over ({circumflex over (r)}₂₂−{circumflex over ({circumflex over (r)}₁₁ mod 2^(N))k₂₁ mod n; and Q=u₂₁*Ĝ₂₂+G.
 20. The method of claim 1 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random number.
 21. The method of claim 3 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random number.
 22. The method of claim 21 wherein the step of deriving the shared secret computes points {circumflex over (Q)}_(B) ₁ =k_(1(A)) ⁻¹Q_(B); {circumflex over (Q)}_(B) ₂ ={circumflex over (r)}_(A){circumflex over (Q)}_(B) ₁ ; and {circumflex over (Q)}_(B) ₃ =(k_(1(A))−k_(2(A))){circumflex over (Q)}_(B) ₁ and then P={circumflex over (Q)}_(B) ₂ +{circumflex over (Q)}_(B) ₃ where the shared secret is Z=x_(P) where x_(P) is the x-coordinate of P, party A uses its transformed random number and party B's public key.
 23. The method of claim 5 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random number.
 24. The method of claim 23 wherein the step of deriving the shared secret computes points {circumflex over (Q)}_(B) ₁ =k_(1(A)) ⁻¹Q_(B); and {circumflex over (Q)}_(B) ₂ ={circumflex over (r)}_(A){circumflex over (Q)}_(B) ₁ and then P={circumflex over (Q)}_(B) ₂ +Q where the shared secret is Z=x_(P) where x_(P) is the x-coordinate of P, party A uses its transformed random number and party B's public key.
 25. The method of claim 11 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random numbers.
 26. The method of claim 25 wherein the step of deriving the shared secret computes {circumflex over (Q)}_(B) ₁₁ =k_(11(A)) ⁻¹Q_(B); {circumflex over (Q)}_(B) ₁₂ =(2{circumflex over ({circumflex over (r)}_(11(A))−{circumflex over ({circumflex over (r)}_(22(A)) mod 2^(N)){circumflex over (Q)}_(B) ₁₁ ; and {circumflex over (Q)}_(B) ₁₃ =(k_(11(A))−k_(12(A))){circumflex over (Q)}_(B) ₁₁ and then P={circumflex over (Q)}_(B) ₁₂ +{circumflex over (Q)}_(B) ₁₃ where the shared secret is Z=x_(P) where x_(P) is the x-coordinate of P, party A uses its transformed random numbers and party B's public key.
 27. The method of claim 13 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random numbers.
 28. The method of claim 27 wherein the step of deriving the shared secret computes {circumflex over (Q)}_(B) ₂₁ =k_(21(A)) ⁻¹Q_(B); {circumflex over (Q)}_(B) ₂₂ ({circumflex over ({circumflex over (r)}−_(22(A))−{circumflex over ({circumflex over (r)}_(11(A)) mod 2^(N)){circumflex over (Q)}_(B) ₂₁ ; and {circumflex over (Q)}_(B) ₂₃ =(k_(21(A))−k_(22(A))){circumflex over (Q)}_(B) ₂₁ and then P={circumflex over (Q)}_(B) ₂₂ +{circumflex over (Q)}_(B) ₂₃ where the shared secret is Z=x_(P) where x_(P) is the x-coordinate of P, party A uses its transformed random numbers and party B's public key.
 29. The method of claim 17 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random number.
 30. The method of claim 29 wherein the step of deriving the shared secret computes {circumflex over (Q)}_(B) ₁₁ =k_(11(A)) ⁻¹Q_(B); and {circumflex over (Q)}_(B) ₁₂ =(2{circumflex over ({circumflex over (r)}_(11(A))−{circumflex over ({circumflex over (r)}_(22(A))mod 2^(N)){circumflex over (Q)}_(B) ₁₁ ; and then P={circumflex over (Q)}_(B) ₁₂ +Q_(B) where the shared secret is Z=x_(P) where x_(P) is the x-coordinate of P, party A uses its transformed random number and party B's public key.
 31. The method of claim 19 further comprising the steps of receiving a public key from another party and deriving a shared secret therefrom using the transformed random number.
 32. The method of claim 31 wherein the step of deriving the shared secret computes {circumflex over (Q)}_(B) ₂₁ =k_(21(A)) ⁻¹Q_(B); and {circumflex over (Q)}_(B) ₂₂ =({circumflex over ({circumflex over (r)}_(22(A))−{circumflex over ({circumflex over (r)}_(11(A)) mod 2^(N)){circumflex over (Q)}_(B) ₂₁ ; and then P={circumflex over (Q)}_(B) ₂₂ +Q_(B) where the shared secret is Z=x_(P) where x_(P) is the x-coordinate of P, party A uses its transformed random numbers and party B's public key.
 33. The method of any claim 1, wherein some values are computed off-line. 